% TEMPLATE for Usenix papers, specifically to meet requirements of
%  USENIX '05
% originally a template for producing IEEE-format articles using LaTeX.
%   written by Matthew Ward, CS Department, Worcester Polytechnic Institute.
% adapted by David Beazley for his excellent SWIG paper in Proceedings,
%   Tcl 96
% turned into a smartass generic template by De Clarke, with thanks to
%   both the above pioneers
% use at your own risk.  Complaints to /dev/null.
% make it two column with no page numbering, default is 10 point

% Munged by Fred Douglis <douglis@research.att.com> 10/97 to separate
% the .sty file from the LaTeX source template, so that people can
% more easily include the .sty file into an existing document.  Also
% changed to more closely follow the style guidelines as represented
% by the Word sample file. 
% This version uses the latex2e styles, not the very ancient 2.09 stuff.
\documentclass[letterpaper,twocolumn,10pt]{article}
\usepackage{usenix,epsfig,endnotes,xspace,color}

% Name candidates:
%  Anza
%  Void 
%  Station (from Genesis's Grand Central component) 
%  TARDIS: Atomic, Recoverable, Datamodel Independent Storage
% EAB: flex, basis, stable, dura
% Stasys:  SYStem for Adaptable Transactional Storage: 

\newcommand{\yad}{Stasis\xspace}
\newcommand{\yads}{Stasis'\xspace}
\newcommand{\oasys}{Oasys\xspace}

\newcommand{\diff}[1]{\textcolor{blue}{\bf #1}}
\newcommand{\eab}[1]{\textcolor{red}{\bf EAB: #1}}
\newcommand{\rcs}[1]{\textcolor{green}{\bf RCS: #1}}
%\newcommand{\mjd}[1]{\textcolor{blue}{\bf MJD: #1}}

\newcommand{\eat}[1]{}

\begin{document}

%don't want date printed
\date{}


%make title bold and 14 pt font (Latex default is non-bold, 16 pt)
\title{\Large \bf \yad: System for Adaptable, Transactional Storage}

%for single author (just remove % characters)
\author{
{\rm Russell Sears}\\
UC Berkeley
\and
{\rm Eric Brewer}\\
UC Berkeley
} % end author

\maketitle

% Use the following at camera-ready time to suppress page numbers.
% Comment it out when you first submit the paper for review.
%\thispagestyle{empty}


%\subsection*{Abstract}

{\em An increasing range of applications require robust support for atomic, durable and concurrent
transactions.  Databases provide the default solution, but force
applications to interact via SQL and to forfeit control over data
layout and access mechanisms.  We argue there is a gap between DBMSs and file systems that limits designers of data-oriented applications.

\yad is a storage framework that incorporates ideas from traditional
write-ahead-logging storage algorithms and file systems.
It provides applications with flexible control over data structures, data layout, performance and robustness properties.
\yad enables the development of
unforeseen variants on transactional storage by generalizing
write-ahead-logging algorithms.  Our partial implementation of these
ideas already provides specialized (and cleaner) semantics to applications.

We evaluate the performance of a traditional transactional storage
system based on \yad, and show that it performs favorably relative to existing
systems.  We present examples that make use of custom access methods, modified
buffer manager semantics, direct log file manipulation, and LSN-free
pages.  These examples facilitate sophisticated performance 
optimizations such as zero-copy I/O.  These extensions are composable,
easy to implement and significantly improve performance.

}
%We argue that our ability to support such a diverse range of
%transactional systems stems directly from our rejection of
%assumptions made by early database designers.  These assumptions
%permeate ``database toolkit'' research.  We attribute the success of
%low-level transaction processing libraries (such as Berkeley DB) to
%a partial break from traditional database dogma.

% entries, and 
% to reduce memory and
%CPU overhead, reorder log entries for increased efficiency, and do
%away with per-page LSNs in order to perform zero-copy transactional
%I/O.  
%We argue that encapsulation allows applications to compose
%extensions.

%These ideas have been partially implemented, and initial performance
%figures, and experience using the library compare favorably with
%existing systems.



\section{Introduction}

As our reliance on computing infrastructure increases, a wider range of 
applications require robust data management.  Traditionally, data management
has been the province of database management systems (DBMSs), which are
well-suited to enterprise applications, but lead to poor support for
systems such as web services, search engines, version systems, work-flow 
applications, bioinformatics, grid computing and scientific computing.  These 
applications have complex transactional storage requirements
but do not fit well
onto SQL or the monolithic approach of current databases.  

Simply providing
access to a database system's internal storage module is an improvement.
However, many of these applications require special transactional properties 
that general purpose transactional storage systems do not provide.  In
fact, DBMSs are often not used for these systems, which instead
implement custom, ad-hoc data management tools on top of file
systems.

A typical example of this mismatch is in the support for
persistent objects.
% in Java, called {\em Enterprise Java Beans}
%(EJB). 
In a typical usage, an array of objects is made persistent by
mapping each object to a row in a table (or sometimes multiple
tables)~\cite{hibernate} and then issuing queries to keep the objects and
rows consistent. An update must confirm it has the current
version, modify the object, write out a serialized version using the
SQL update command and commit.  Also, for efficiency, most systems must 
buffer two copies of the application's working set in memory.  
This is an awkward and slow mechanism.

Bioinformatics systems perform complex scientific
computations over large, semi-structured databases with rapidly evolving schemas.  Versioning and
lineage tracking are also key concerns.  Relational databases support
none of these requirements well.  Instead, office suites, ad-hoc
text-based formats and Perl scripts are used for data management~\cite{perl} (with mixed success~\cite{excel}).

\eat{
Examples of real world systems that currently fall into this category
are web search engines, document repositories, large-scale web-email
services, map and trip planning services, ticket reservation systems,
photo and video repositories, bioinformatics, version control systems,
work-flow applications, CAD/VLSI applications and directory services.

In short, we believe that a fundamental architectural shift in
transactional storage is necessary before general purpose storage
systems are of practical use to modern applications.
Until this change occurs, databases' imposition of unwanted
abstraction upon their users will restrict system designs and
implementations.
}

%In short, reliable data management has become as unavoidable as any
%other operating system service.  As this has happened, database
%designs have not incorporated this decade-old lesson from operating
%systems research:
%
%\begin{quote} The defining tragedy of the operating systems community
%  has been the definition of an operating system as software that both
%  multiplexes and {\em abstracts} physical resources...The solution we
%  propose is simple: complete elimination of operating systems
%  abstractions by lowering the operating system interface to the
%  hardware level~\cite{engler95}.
%\end{quote}

%The widespread success of lower-level transactional storage libraries
%(such as Berkeley DB) is a sign of these trends.  However, the level
%of abstraction provided by these systems is well above the hardware
%level, and applications that resort to ad-hoc storage mechanisms are
%still common.

This paper presents \yad, a library that provides transactional
storage at a level of abstraction as close to the hardware as
possible.  The library can support special purpose, transactional
storage interfaces in addition to ACID database-style interfaces to
abstract data models.  \yad incorporates techniques from databases
(e.g. write-ahead-logging) and systems (e.g. zero-copy techniques).
Our goal is to combine the flexibility and layering of low-level
abstractions typical for systems work with the complete semantics
that exemplify the database field.

By {\em flexible} we mean that \yad{}  can implement a wide
range of transactional data structures, that it can support a variety
of policies for locking, commit, clusters and buffer management.
Also, it is extensible for new core operations
and new data structures. It is this flexibility that allows the
support of a wide range of systems.

By {\em complete} we mean full redo/undo logging that supports
both {\em no force}, which provides durability with only log writes,
and {\em steal}, which allows dirty pages to be written out prematurely
to reduce memory pressure. By complete, we also
mean support for media recovery, which is the ability to roll
forward from an archived copy, and support for error-handling,
clusters, and multithreading. These requirements are difficult
to meet and form the {\em raison d'\^etre} for \yad{}: the framework
delivers these properties as reusable building blocks for systems
that implement complete transactions.

Through examples and their good performance, we show how \yad{}
supports a wide range of uses that fall in the gap between 
database and filesystem technologies, including
persistent objects, graph or XML based applications, and recoverable
virtual memory~\cite{lrvm}.  

For example, on an object serialization workload, we provide up to 
a 4x speedup over an in-process 
MySQL implementation and a 3x speedup over Berkeley DB while 
cutting memory usage in half (Section~\ref{sec:oasys}). 

We implemented this extension in 150 lines of C, including comments and boilerplate.  We did not have this type of optimization
in mind when we wrote \yad.  In fact, the idea came from a potential 
user that is not familiar with \yad.

%\e ab{others?  CVS, windows registry, berk DB, Grid FS?}
%\r cs{maybe in related work?}

This paper begins by contrasting \yads approach with that of
conventional database and transactional storage systems.  It proceeds
to discuss write-ahead-logging, and describe ways in which \yad can be
customized to implement many existing (and some new) write-ahead-logging variants.  Implementations of some of these variants are
presented, and benchmarked against popular real-world systems.  We
conclude with a survey of the technologies the \yad implementation is
based upon.

An (early) open-source implementation of
the ideas presented here is available.

\section{\yad is not a Database}
\label{sec:notDB}
Database research has a long history, including the development of
many technologies that our system builds upon.  This section explains
why databases are fundamentally inappropriate tools for system
developers.  The problems we present here have been the focus of
database systems and research projects for at least 25 years.

\subsection{The database abstraction}

Database systems are often thought of in terms of the high-level
abstractions they present.  For instance, relational database systems
implement the relational model~\cite{codd}, object oriented
databases implement object abstractions, XML databases implement
hierarchical datasets, and so on.  Before the relational model,
navigational databases implemented pointer- and record-based data models.

An early survey of database implementations sought to enumerate the
fundamental components used by database system implementors.  This
survey was performed due to difficulties in extending database systems
into new application domains.  It divided internal database
routines into two broad modules: {\em conceptual
mappings}~\cite{batoryConceptual} and {\em physical
database models}~\cite{batoryPhysical}.

%A physical model would then translate a set of tuples into an
%on-disk B-Tree, and provide support for iterators and range-based query
%operations.

It is the responsibility of a database implementor to choose a set of
conceptual mappings that implement the desired higher-level
abstraction (such as the relational model).  The physical data model
is chosen to efficiently support the set of mappings that are built on
top of it.

\diff{A conceptual mapping based on the relational model might
translate a relation into a set of keyed tuples.  If the database were
going to be used for short, write-intensive and high-concurrency
transactions (OLTP), the physical model would probably translate sets
of tuples into an on-disk B-Tree.  In contrast, if the database needed
to support long-running, read only aggregation queries (OLAP), a
physical model tuned for such queries\rcs{be more concrete here} would
be more appropriate.  While both OLTP and OLAP databases are based
upon the relational model they make use of different physical models
in order to serve different classes of applications.}

A key observation of this paper is that no known physical data model
can efficiently support more than a small percentage of today's applications.  

Instead of attempting to create such a model after decades of database
research has failed to produce one, we opt to provide a transactional
storage model that mimics the primitives provided by modern hardware.
This makes it easy for system designers to implement most of the data
models that the underlying hardware can support, or to
abandon the database approach entirely, and forgo the use of a
structured physical model or abstract conceptual mappings.

\subsection{Extensible transaction systems} 
\label{sec:otherDBs}
This section contains discussion of database systems with goals similar to ours.
Although these projects were
successful in many respects, they fundamentally aimed to implement a
extensible data model, rather than build transactions from the bottom up.
In each case, this limits the applicability of their implementations.

\subsubsection{Extensible databases}

Genesis~\cite{genesis}, an early database toolkit, was built in terms
of a physical data model and the conceptual mappings described above.
It is designed to allow database implementors to easily swap out
implementations of the various components defined by its framework.
Like subsequent systems (including \yad), it allows its users to
implement custom operations.

Subsequent extensible database work builds upon these foundations.
The Exodus~\cite{exodus} database toolkit is the successor to
Genesis. It supports the automatic generation of query optimizers and
execution engines based upon abstract data type definitions, access
methods and cost models provided by its users.

Although further discussion is beyond the scope of this paper,
object-oriented database systems and relational databases with
support for user-definable abstract data types (such as in
Postgres~\cite{postgres}) were the primary competitors to extensible
database toolkits.  Ideas from all of these systems have been
incorporated into the mechanisms that support user-definable types in
current database systems.

One can characterize the difference between database toolkits and
extensible database servers in terms of early and late binding.  With
a database toolkit, new types are defined when the database server is
compiled.  In today's object-relational database systems, new types
are defined at runtime.  Each approach has its advantages.  However,
both types of systems aim to extend a high-level data model with new 
abstract data types, and thus are quite limited in the range of new 
applications they support.  In hindsight, it is not surprising that this kind of 
extensibility has had little impact on the range of applications 
we listed above.

\subsubsection{Berkeley DB}

%System R was one of the first relational database implementations, and
%defined a clean separation between its query processor and its storage
%subsystem.  In fact, it supported a simple navigational interface to
%the storage subsystem, which remains the architecture for modern
%databases.

Berkeley DB is a highly successful alternative to conventional
databases.  At its core, it provides the physical database
(relational storage system) of a conventional database server.
%It is based on the
%observation that the storage subsystem is a more general (and less
%abstract) component than a monolithic database, and provides a
%stand-alone implementation of the storage primitives built into 
%most relational database systems~\cite{libtp}.  
In particular, 
it provides fully transactional (ACID) operations over B-Trees, 
hashtables, and other access methods.  It provides flags that 
let its users tweak various aspects of the performance of these
primitives, and selectively disable the features it provides~\cite{libtp}.

With the
exception of the benchmark designed to fairly compare the two systems, none of the \yad 
applications presented in Section~\ref{sec:extensions} are efficiently
supported by Berkeley DB.   This is a result of Berkeley DB's  
assumptions regarding workloads and decisions regarding low level data
representation.  Thus, although Berkeley DB could be built on top of \yad,
Berkeley DB's data model, and write-ahead-logging system are too specialized to support \yad.



%cover P2 (the old one, not Pier 2 if there is time...

\subsubsection{Better databases}

The database community is also aware of this gap. 
A recent survey~\cite{riscDB} enumerates problems that plague users of
state-of-the-art database systems, and finds that database implementations fail to support the
needs of modern applications.  Essentially, it argues that modern 
databases are too complex to be implemented (or understood) 
as a monolithic entity.

It supports this argument with real-world evidence that suggests
database servers are too unpredictable and unmanagable to
scale up the size of today's systems.  Similarly, they are a poor fit
for small devices.  SQL's declarative interface only complicates the
situation.

%In large systems, this manifests itself as
%manageability and tuning issues that prevent databases from predictably
%servicing diverse, large scale, declarative, workloads.  
%On small devices, footprint, predictable performance, and power consumption are
%primary concerns that database systems do not address.

%The survey argues that these problems cannot be adequately addressed without a fundamental shift in the architectures that underly database systems.  Complete, modern database
%implementations are generally incomprehensible and
%irreproducible, hindering further research.  
The study concludes 
by suggesting the adoption of {\em RISC} database architectures, both as a resource for researchers and as a 
real-world database system.

RISC databases have many elements in common with
database toolkits.  However, they take the database toolkit idea one
step further, and suggest standardizing the interfaces of the
toolkit's internal components, allowing multiple organizations to
compete to improve each module.  The idea is to produce a research
platform that enables specialization and shares the effort required to build a full database~\cite{riscDB}.

We agree with the motivations behind RISC databases, and to build 
databases from interchangeable modules exists.  In fact, is our hope
that our system will mature to the point where it can support 
a competitive relational database.  However this is
not our primary goal.  
%Instead, we are interested in supporting applications that derive
%little benefit from database abstractions, but that need reliable
%storage.  Therefore, 
Instead of building a modular database, we seek
to build a system that enables a wider range of data management options.

%For example, large scale application such as web search, map services,
%e-mail use databases to store unstructured binary data, if at all.

%More recently, WinFS, Microsoft's database based
%file meta data management system, has been replaced in favor of an
%embedded indexing engine that imposes less structure (and provides
%fewer consistency guarantees) than the original
%proposal~\cite{needtocitesomething}.

%Scaling to the very large doesn't work (SAP used DB2 as a hash table
%for years), search engines, cad/VLSI didn't happen.  scalable GIS
%systems use shredded blobs (terraserver, google maps), scaling to many
%was more difficult than implementing from scratch (winfs), scaling
%down doesn't work (variance in performance, footprint),

\section{Conventional Transactions in \yad}

\rcs{This whole section is new, and is intended to replace what is now section 4.}

\rcs{This section is missing references to prior work.  Bill mentioned
PhD theses that talk about this layering, but I've been too busy
coding to read them.}

This section describes how \yad implements transactions that are
similar to those provided by relational database systems.  In addition
to providing a review of how modern transactional systems function,
this section lays out the functionality that \yad provides to the
operations built on top of it.  It also explains how \yads
operations are roughly structured as two levels of abstraction.  

The lower level of a \yad operation provides atomic
updates to regions of the disk.  These updates do not have to deal
with concurrency, but the portion of the page file that they read and
write must be atomically updated, even if the system crashes.

The higher level atomically applies operations
to the page file to provide operations that span multiple pages and
copes with concurrency issues.  Surprisingly, the implementations
of these two layers are only loosely coupled.

Finally, this section describes how \yad manages transaction-duration
locks and discusses the alternatives \yad provides to application developers.

\subsection{Atomic page file operations}

Transactional storage algorithms work because they are able to
atomically update portions of durable storage.  These small atomic
updates are used to bootstrap transactions that are too large to be
applied atomically.  In particular, write ahead logging (and therefore
\yad) relies on the ability to atomically write entries to the log
file.

\subsubsection{Hard drive behavior during a crash}
In practice, a write to a disk page is not atomic.  Two common failure
modes exist.  The first occurs when the disk writes a partial sector
to disk during a crash.  In this case, the drive maintains an internal
checksum, detects a mismatch, and reports it when the page is read.
The second case occurs because pages span multiple sectors.  Drives
may reorder writes on sector boundaries, causing an arbitrary subset
of a page's sectors to be updated during a crash.  

{\em Torn page detection} can be used to detect this phenomonon.  Torn
and corrupted pages may be recovered by restoring the page from
backup.  For simplicity, this section ignores mechanisms that detect
and restore torn pages, and assumes that page writes are atomic.
While the techniques described in this section rely on the ability to
atomically update disk pages, this restriction is relaxed by other 
recovery mechanisms.


\subsubsection{Extending \yad with new operations}

Figure~\ref{fig:structure} shows how custom operations interact with
\yad.  If an application does not need to make use of concurrent
transactions, directly manipulating the page file is as simple as
ensuring that each update to the page file occurs inside of an
operation's implementation.  Operation implementations must be invoked
by registering a callback with \yad at startup, and then calling {\em
Tupdate()} to invoke the operation at runtime.  Each operation should
be deterministic, provide an inverse, and acquire all of its arguments
from a struct that is passed via Tupdate().  (Operations that affect
more than one page, and ones that do not provide inverses will be
described later.)  The same callbacks are used during forward opertion
as during recovery.  Therefore operations provide a single redo
function and a single undo function.  (There is no ``do''
function.)  This reduces the amount of recovery-specific code in the
system.

\subsubsection{\yads Recovery Algorithm}

Recovery relies upon the fact that each log entry is assigned a {\em
Log Sequence Number (LSN)}.  The LSN is monitonically increasing and
unique.  The LSN of the log entry that was most recently applied to
each page is stored with the page, allowing recovery to selectively
replay log entries.  This only works if log entries change exactly one
page, and if they are applied to the page atomically.

Recovery occurs in three phases, Analysis, Redo and Undo.
``Analysis'' is beyond the scope of this paper.  ``Redo'' plays the
log forward in time, applying any updates that did not make it to disk
before the system crashed.  ``Undo'' runs the log backwards in time,
only applying portions that correspond to aborted transactions.  This
section only considers physical undo.  Section~\ref{sec:nta} describes
the distinction between physical and logical undo, and describes
logical undo.  A summary of the stages of recovery and the invariants
they establish is presented in Figure~\ref{fig:conventional-recovery}.

Redo is the only phase that makes use of LSN's stored on pages.
It simply compares the page LSN to the LSN of each log entry.  If the
log entry's LSN is higher than the page LSN, then the log entry is
applied.  Otherwise, the log entry is skipped.  Redo does not write
log entries to disk, as it is replaying events that have already been
recorded.  

However, Undo does write log entries.  In order to prevent repeated
crashes during recovery from causing the log to grow excessively, the
entries that Undo writes tell future invocations of Undo to skip
portions of the transaction that have already been undone.  These log
entries are usually called {\em Compensation Log Records (CLR's)}.
Note that CLR's only cause Undo to skip log entries.  Redo will apply
log entries protected by the CLR, guaranteeing that those updates are
applied to the page file.

There are many other schemes for page level recovery that we could
have chosen.  The scheme desribed above has two particularly nice
properties.  First, pages that were modified by active transactions
may be {\em stolen}; they may be written to disk before a transaction
completes.  This allows transactions to use more memory than is
physically available, and makes it easier to flush frequently written
pages to disk.  Second, pages do not need to be {\em forced}; a
transaction commits simply by flushing the log.  If it had to force
pages to disk it would incur the cost of random I/O.  Also, if
multiple transactions commit in a small window of time, the log only
needs to be forced to disk once.

\subsubsection{Alternatives to Steal / no-Force}

Note that the Redo phase of recovery allows \yad to avoid forcing
pages to disk, while Undo allows pages to be stolen.  For some
applications, the overhead of logging information for Redo or Undo may
outweigh their benefits.  \yads logging discipline provides a simple
solution to this problem.  If a special-purpose operation wants to
avoid writing either the Redo or the Undo information to the log then
it can have the buffer manager pin the page or flush it at commit, and
simply omit the pertinent information from the log entries it
generates.

Recovery's Undo and Redo phases both will process the log entry, but
one of them will have no effect.  If an operation chooses not to
provide a Redo implementation, then its Undo implementation will need
to determine whether or not the Redo was applied.  If it omits Undo,
then Redo must consult recovery to see if it is part of a transaction that
committed.

\subsection{Concurrent Transactions}

Two factors make it more difficult to write operations that may be
used in concurrent transactions.  The first is familiar to anyone that
has written multi-threaded code: Accesses to shared data structures
must be protected by latches (mutexes).  The second problem stems from
the fact that concurrent transactions prevent abort from simply
rolling back the physical updates that a transaction made.
Fortunately, it is straightforward to reduce this second,
transaction-specific, problem to the familiar problem of writing
multi-threaded software.

To understand why redo cannot simply revert each page to the state it
was in before a transaction began, consider an operation that inserts
data into a tree, and the following sequence of events: 
\begin{itemize}
  \item Transaction A inserts data, causing a node to be
    split
  \item Transaction B inserts data into one of the newly
    created nodes
  \item Transaction A calls abort
\end{itemize}
If abort simply restored the pages to the state they were in before A
updated them, then the data item that transaction B inserted would be
lost.  Operations that apply changes to pages without an understanding
of the data they manipulate are called {\em physical operations}.

If we constrain the tree structure to fit on a single page, then the
``insert'' operation's inverse could be a ``remove'' operation.  Such
operations are called {\em logical operations}.  In this case, both
operations would traverse tree nodes to determine what updates should
be applied and modify the tree accordingly.  This would allow abort to
remove A's data from the tree without losing B's updates.  

The problem becomes more complex if we allow the tree to span multiple
pages.  If we use a single log entry to record the update and the
system crashes, then there is no guarantee that the LSNs of the pages
that the log entry manipulates will match, or that the two pages will
contain physically consistent portions of the tree structure.

In general, physical operations cause concurrent transactions to
violate the physical consistency of data structures during abort.
Logical operations that span more than one page cannot safely be
redone during recovery.

{\em Nested Top Actions} provide an elegant solution to this problem.
A nested top action uses physical undo while a data structure is being
upated, and then atomically switches to logical undo once the data
structure is internally consistent.  Nested top actions work by
performing physical operations on a data structure, and then
registering a CLR.  The CLR contains a logical undo entry for the
operation.  When recovery and abort encounter a CLR they skip the
physical undo entries, and instead apply the logical undo.

From the perspective of an operation implementation, a nested top
action protects logical undo functions from seeing temporary
inconsistencies introduced by operations that span pages.  Since
latches protect other threads from the same set of inconsistencies,
the proper use of nested top actions is similar to the development of
thread-safe code.

This leads to a mechanical approach that converts non-reentrant
operations that do not support concurrent transactions into reentrant,
concurrent operations:

\begin{enumerate}
\item Wrap a mutex around each operation.  With care, it may be possible to use finer-grained locks, but it is rarely necessary.
\item Define a {\em logical} UNDO for each operation (rather than just
  using a set of page-level UNDO's).  For example, this is easy for a
  hashtable: the UNDO for {\em insert} is {\em remove}.  \diff{This logical
  undo function should arrange to acquire the mutex when invoked by
  abort or recovery.}
\item Add a ``begin nested
  top action'' right after the mutex acquisition, and an ``end
  nested top action'' right before the mutex is released.  \diff{\yad provides a default nested top action implementation as an extension.}
\end{enumerate}

If the transaction that encloses a nested top action aborts, the
logical undo will {\em compensate} for the effects of the operation,
leaving structural changes intact.  If a transaction should perform
some action regardless of whether or not it commits, a nested top
action with a ``no-op'' as its inverse is a convenient way of applying
the change.  Nested top actions do not cause the log to be forced to disk, so
such changes will not be durable until the log is manually forced, or
until the updates eventually reach disk.

This section described how concurrent, thread-safe operations can be
developed.  These operations provide building blocks for concurrent
transactions, and are fairly easy to develop.  Interestingly, any
mechanism that applies atomic physical updates to the page file can be
used as the basis of a nested top action.  However, concurrent
operations are of little help if an application is not able to safely
combine them to create concurrent transactions.  

\subsection{Application-specific Locking}

Note that the transactions described above only provide the
``Atomicity'' and ``Durability'' properties of ACID.  ``Isolation'' is
typically provided by locking, which is a higher-level (but
comaptible) layer.  ``Consistency'' is less well defined but comes in
part from low-level mutexes that avoid races, and partially from
higher level constructs such as unique key requirements.  \yad
supports this by distinguishing between {\em latches} and {\em locks}.
Latches are provided using operating system mutexes, and are held for
short periods of time.  \yads default data structures use latches in a
way that avoids deadlock.  This section will describe the latching
protocols that \yad makes use of, and describes two custom lock
managers that \yads allocation routines use to implement layout policies and provide deadlock avoidance.

This allows higher level code to treat \yad as a conventional
reentrant data structure library.  It is the application's
responsibility to provide locking, whether it be via a database-style
lock manager, or an application-specific locking protocol.  Note that
locking schemes may be layered.  For example, when \yad allocates a
record, it first calls a region allocator that allocates contiguous
sets of pages, and then it allocates a record on one of those pages.

The record allocator and the region allocator each contain custom lock
management.  If transaction A frees some storage, transaction B reuses
the storage and commits, and then transaction A aborts, then the
storage would be double allocated.  The region allocator (which is
infrequently called, and not concerned with locality) records the id
of the transaction that created a region of freespace, and does not
coalesce or reuse any storage associated with an active transaction.

On the other hand, the record allocator is called frequently, and is
concerned with locality.  Therefore, it associates a set of pages with
each transaction, and keeps track of deallocation events, making sure
that space on a page is never over reserved.  Providing each
transaction with a seperate pool of freespace should increase
concurrency and locality.  This allocation strategy was inspired by
Hoard, a malloc implementation for SMP machines.

Note that both lock managers have implementations that are tied to the
code they service, both implement deadlock avoidance, and both are
transparent to higher layers.  General purpose database lock managers
provide none of these features, supporting the idea that special
purpose lock managers are a useful abstraction.\rcs{This would be a
good place to cite Bill and others on higher level locking protocols}

Locking is largely orthoganol to the concepts desribed in this paper.
We make no assumptions regarding lock managers being used by higher
level code in the remainder of this discussion.

\section{LSN-free pages.}
\label{sec:lsn-free}
\rcs{After working through the torn page argument, I realized that
this style of transaction allows you to produce log entries that make
non-localized updates to the page file.  I think this means that we
can avoid writing out physical undo information for our nested top
actions, and write the redo in a single entry.  This is particularly
interesting because LSN-free recovery will break horribly if the
logical undo of a nested transaction reads or writes bytes that happen
to be written or read by a partial nested top action that is being
physically rolled back.  If nested top actions are atomic log
entities, the problem cannot occur.  Of course, this approach has its
limits; the longer a log entry is, the more transactions will block
waiting for it to be appended to the end of the log.  Also, if we mix
LSN and LSN-free operations in the same nested top action, we need the
physical undo.}

\rcs{ Do something with this text... Logical operations that are constrained to a single page are often
called {\em physiological operations}, and are used throughout \yad.
Note that physioloical operations are not necessarily idempotent, and
they rely upon the consistency of the page they modify.  In
Section~\ref{XXX}, \yad used page LSN's to guarantee that the
operations recorded in the log are atomically applied exactly
once. The recovery scheme described in this section does not provide
these guarantees and is incompatible with physiological operations.}

The recovery algorithm described above uses LSN's to determine the
version number of each page during recovery.  This is a common
technique.  As far as we know, is used by all database systems that
update data in place.  Unfortunately, this makes it difficult to map
large objects onto pages, as the LSN's break up the object.  It
is tempting to store the LSN's elsewhere, but then they would not be
written atomically with their page, which defeats their purpose.  

This section explains how we can avoid storing LSN's on pages in \yad
without giving up durable transactional updates.  In the process, we
are able to relax the atomicity assumptions that we make regarding
writes to disk.  These relaxed assumptions allow recovery to repair
torn pages without performing media recovery, and allow arbitrary
ranges of the page file to be updated by a single physical operation.

\yads implementation does not currently support the recovery algorithm
described in this section.  However, \yad avoids hard-coding most of
the relevant subsytems.  LSN-free pages are essentially an alternative
protocol for atomically and durably applying updates to the page file.
We plan to eventually support the coexistance of LSN-free pages,
traditional pages, and similar third-party modules within the same
page file, log, transactions, and even logical operations.

\subsection{Blind writes}
Recall that LSN's were introduced to prevent recovery from applying
updates more than once, and to prevent recovery from applying old
updates to newer versions of pages.  This was necessary because some
operations that manipulate pages are not idempotent, or simply make
use of state stored in the page.  We can avoid such problems by
eliminating such operations and instead making use of deterministic
REDO operations that do not examine page state.  We call such
operations ``blind writes.''

For concreteness, assume that all physical operations produce log
entries that contain a set of byte ranges, and the pre- and
post-value of each byte in the range.  

Recovery works the same way as it does above, except that is computes
a lower bound of each page LSN instead of reading the LSN from the
page.  One possible lower bound is the LSN of the most recent log
truncation or checkpoint.  Alternatively, \yad could occasionally
write information about the state of the buffer manager to the log.

Although the mechanism used for recovery is similar, the invariants
maintained during recovery have changed.  With conventional
transactions, if a page in the page file is internally consistent
immediately after a crash, then the page will remain internally
consistent throughout the recovery process.  This is not the case with
our LSN-free scheme.  If a consistent, relatively new, version of a
page is on disk immediately after a crash, then that page may be
overwritten using physical log entries that are older than it.
Therefore, the page will contain a mixture of new and old bytes, and
any data structures stored on the page may be inconsistent.  However,
once the redo phase is complete, any old bytes will be overwritten by
their most recent values, so the page will contain an internally
consistent, up-to-date version of itself.
(Section~\ref{sec:torn-page} explains this in more detail.)

Undo can then proceed normally as long as the operations that it logs
to disk only perform blind-writes.  Since this restriction also
applies to normal operations, we suspect this will not pose many
practical problems.

As long as operations are limited to blind writes we do not need to
store LSN's in pages.  The rest of this section describes how this
allows standard filesystem and database optimizations to be easily
combined, and shows that the removal of LSN's from pages actually
simplifies some aspects of recovery.

\subsection{Zero-copy I/O} 

We originally developed LSN-free pages as an efficient method for
storing large (multi-page) objects in the filesystem.  If a large
object is stored in pages that contain LSN's, then in order to read
that large object the system must read each page individually, and
then use the CPU to copy the portions of the page that contain data
into a second buffer.  

Compare
this approach to a modern filesystem, which allows applications to
perform a DMA copy of the data into memory, avoiding the expensive
byte-by-byte copy of the data, and allowing the CPU to be used for
more productive purposes.  Furthermore, modern operating systems allow
network services to use DMA and network adaptor hardware to read data
from disk, and send it over a network socket without passing it
through the CPU.  Again, this frees the CPU, allowing it to perform
other tasks.

We believe that LSN free pages will allow reads to make use of such
optimizations in a straightforward fashion.  Zero copy writes are more challenging, but could be
performed by performing a DMA write to a portion of the log file.
However, doing this complicates log truncation, and does not address
the problem of updating the page file.  We suspect that contributions
from the log based filesystem~\cite{lfs} literature can address these problems in
a straightforward fashion.  In particular, we imagine storing 
portions of the log (the portion that stores the blob) in the 
page file, or other addressable storage.  In the worst case, 
the blob would have to be relocated in order to defragment the 
storage.  Assuming the blob was relocated once, this would amount 
to a total of three, mostly sequential disk operations.  (Two 
writes and one read.)  However, in the best case, the blob would only need to written once.
In contrast, a conventional atomic blob implementation would always need 
to write the blob twice.

Alternatively, we could use DMA to overwrite the blob in the page file
in a non-atomic fashion, providing filesystem style semantics.
(Existing database servers often provide this mode based on the
observation that many blobs are static data that does not really need
to be updated transactionally.~\cite{sqlserver}) Of course, \yad could
also support other approaches to blob storage, such as B-Tree layouts
that allow arbitrary insertions and deletions in the middle of
objects~\cite{esm}.

\subsection{Concurrent recoverable virtual memory}

Our LSN-free pages are somewhat similar to the recovery scheme used by
RVM, recoverable virtual memory.  That system used purely physical
logging and LSN-free pages so that it could use mmap() to map portions
of the page file into application memory\cite{lrvm}.  However, without
support for logical log entries and nested top actions, it would be
difficult to implement a concurrent, durable data structure using RVM.
We plan to add RVM style transactional memory to \yad in a way that is
compatible with fully concurrent collections such as hash tables and
tree structures.  Of course, since \yad will support coexistance of
conventional and LSN-free pages, applications would be free to use the
\yad data structure implementations as well.

\subsection{Page-independent transactions}
\rcs{I don't like this section heading...}  Recovery schemes that make
use of per-page LSN's assume that each page is written to disk
atomically even though that is generally not the case.  Such schemes
deal with this problem by using page formats that allow partially
written pages to be detected.  Media recovery allows them to recover
these pages.

The Redo phase of the LSN-free recovery algorithm actually creates a
torn page each time it applies an old log entry to a new page.
However, it guarantees that all such torn pages will be repaired by
the time Redo completes.  In the process, it also repairs any pages
that were torn by a crash.  Instead of relying upon atomic page
updates, LSN-free recovery relies upon a weaker property.

For LSN-free recovery to work properly after a crash, each bit in
persistent storage must be either:

\begin{enumerate}
\item The old version of a bit that was being overwritten during a crash.
\item The newest version of the bit written to storage.
\item Detectably corrupt (the storage hardware issues an error when the
  bit is read).
\end{enumerate}

Modern drives provide these properties at a sector level: Each sector
is atomically updated, or it fails a checksum when read, triggering an
error.  If a sector is found to be corrupt, then media recovery can be
used to restore the sector from the most recent backup.

Figure~\ref{fig:todo} provides an example page, and a number of log
entries that were applied to it.  Assume that the initial version of
the page, with LSN $0$, is on disk, and the disk is in the process of
writing out the version with LSN $2$ when the system crashes.  When
recovery reads the page from disk, it may encounter any combination of
sectors from these two versions.

Note that the first and last two sectors are not overwritten by any
of the log entries that Redo will play back.  Therefore, their value
is unchanged in both versions of the page.  Since Redo will not change
them, we know that they will have the correct value when it completes.
The remainder of the sectors are overwritten at some point in the log.
If we constrain the updates to overwrite an entire page at once, then
the initial on-disk value of these sectors would not have any affect
on the outcome of Redo.  Furthermore, since the redo entries are
played back in order, each sector would contain the most up to date
version after redo.

Of course, we do not want to constrain log entries to update entire
sectors at once.  In order to support finer grained logging, we simply
repeat the above argument on the byte or bit level.  Each bit is
either overwritten by redo, or has a known, correct, value before
redo.  Since all operations performed by redo are blind writes, they
can be applied regardless of whether the page is logically consistent.

Since LSN-free recovery only relies upon atomic updates at the bit
level, it prevents pages from becoming a limit to the size of atomic
page file updates.  This allows operations to atomically manipulate
(potentially non-contiguous) regions of arbitrary size by producing a
single log entry.

This is particularly convenient when dealing with nested top actions.
Normally, a nested top action performs a number of updates to the page
file, and logs a physical undo entry for each one.  Upon completion,
it writes a logical undo entry.  The physical undo entries take up
space in the log, and reduce the amount of log bandwidth available for
other tasks.  In cases where a nested top action can be completed by
only logging blind writes, the logical undo that would normally
complete the nested top action can replace the physical undo entries.
This only works because the log entry and its logical undo are
atomically applied to the page file.  With conventional transactions,
this technique is limited to operations that update a single page.
LSN-free pages remove this limitation.

\section{Transactional Pages}

\rcs{I plan to cut out all of section 4 as it currently exists, but it
still contains stuff that needs to be in section 3.}

\rcs{I think we should avoid the term ``transactional pages''.  In the
LSN-free pages discussion, we rely upon the atomicity of the
application of each log operation after REDO.  We should say that we
will start by talking about updates that are within a single page (and
assumed to be applied to disk atomically), but that there are other
ways to atomically update storage.  Multi-page transactions break the
atomicity assumption because their results are not applied to disk
atomically.  Concurrent transactions break the assumption that a
series of physical undos is the inverse of a transaction.  Nested top
actions restore these two broken invariants, but are orthoganol to the
mechanisms that apply the atomic updates.  (This is why LSN free pages
are compatible with Nested Top Actions.)  I think this section should
do three things.  (1)Explain the distinction between lower level
atomicity and nested top actions, (2) Explain how recovery works. (3)
Tease out aspects of transactional storage that recovery doesn't need
to worry about.}

Section~\ref{sec:notDB} described the ways in which a top-down data model
limits the generality and flexibility of databases.  In this section,
we cover the basic bottom-up approach of \yad: {\em transactional
pages}. Although similar to the underlying write-ahead-logging
approaches of databases, particularly ARIES~\cite{aries}, \yads
bottom-up approach yields unexpected flexibility.

Transactional pages provide the properties of transactions, but
only allow updates within a single page in the simplest case.  After
covering the single-page case, we explore multi-page transactions,
which enable a complete transaction system.

In this model, pages are the in-memory representation of disk blocks
and thus must be the same size.  Pages are a convenient abstraction
because the write back of a page (disk block) is normally atomic,
giving us a foundation for larger atomic actions. In practice, disk
blocks are not always atomic, but the disk can detect partial writes
via checksums.  Thus, we actually depend only on detection of
non-atomicity, which we treat as media failure.  One nice property of
\yad is that we can roll forward an individual page from an archive copy to
recover from media failures.\rcs{Torn page detection...}

A subtlety of transactional pages is that they technically only
provide the ``atomicity'' and ``durability'' of ACID
transactions.\endnote{The ``A'' in ACID really means atomic persistence
of data, rather than atomic in-memory updates, as the term is normally
used in systems work; %~\cite{GR97}; 
the latter is covered by ``C'' and
``I''.}  This is because ``isolation'' comes typically from locking, which
is a higher (but compatible) layer. ``Consistency'' is less well defined
but comes in part from transactional pages (from mutexes to avoid race
conditions), and in part from higher layers (e.g. unique key
requirements). To support these, \yad distinguishes between {\em
latches} and {\em locks}.  A latch corresponds to an OS mutex, and is
held for a short period of time.  All of \yads default data structures
use latches in a way that avoids deadlock. This allows
multithreaded code to treat \yad as a conventional reentrant data structure
library.  Applications that want conventional isolation
(serializability) can make use of a lock manager.

\eat{
\yad uses write-ahead-logging to support the
four properties of transactional storage: Atomicity, Consistency,
Isolation and Durability.  Like existing transactional storage systems,
\yad allows applications to disable or choose different variants of each
property.

However, \yad takes customization of transactional semantics one step
further, allowing applications to add support for transactional
semantics that we have not anticipated.  We do not believe that
we can anticipate every possible variation of write-ahead-logging.  
However, we
have observed that most changes that we are interested in making
involve a few common underlying primitives.  

As we have
implemented new extensions, we have located portions of the system
that are prone to change, and have extended the API accordingly.  Our
goal is to allow applications to implement their own modules to
replace our implementations of each of the major write-ahead-logging
components.
}


\subsection{Single-page Transactions}

In this section we show how to implement single-page transactions.
This is not at all novel, and is in fact based on ARIES~\cite{aries},
but it forms important background.  We also gloss over many important
and well-known optimizations that \yad exploits, such as group
commit.%~\cite{group-commit}.  
These aspects of recovery algorithms are
described in the literature, and in any good textbook that describes
database implementations.  They are not particularly important to our
discussion, so we do not cover them.

The trivial way to achieve single-page transactions is simply to apply
all the updates to the page and then write it out on commit. The page
must be pinned until the transaction commits to avoid ``dirty'' data
(uncommitted data on disk), but no logging is required.  As disk
block writes are atomic, this ensures that we provide the ``A'' and ``D''
of ACID.

This approach scales poorly to multiple pages since we must {\em force} pages to disk
on commit and wait for a (random access) synchronous write to
complete. By using a write-ahead log, we can support {\em no force}
transactions: we write (sequential) ``redo'' information to the log on commit, and
then can write the  pages later. If we crash, we can use the log to
redo the lost updates during recovery.

For this to work, recovery must be able to decide which updates to
re-apply.  This is solved by using a per-page sequence number called a
{\em log sequence number \diff{(LSN)}}. Each log entry contains the sequence
number, and each page contains the sequence number of the last applied
update.  Thus on recovery, we load a page, look at its sequence
number, and re-apply all later updates.  Similarly, to restore a page
from archive we use the same process, but with likely many more
updates to apply.

We also need to make sure that only the results of committed
transactions still exist after recovery.  This is best done by writing
a commit record to the log during the commit.  If the system pins uncommitted
dirty pages in memory, recovery does not need to worry about undoing 
any updates.  Therefore recovery simply plays back unapplied redo records from 
transactions that have commit records.

However, pinning the pages of active transactions in memory is problematic.
First, a single transaction may need more pages than can be pinned at
one time. Second, under concurrent transactions, a given page may be
pinned forever as long as it has at least one active transaction in
progress all the time.  To avoid these problems, transaction systems
support {\em steal}, which means that pages can be written back
before a transaction commits. 

Thus, on recovery a page may contain data that never committed and the
corresponding updates must be rolled back.  To enable this, ``undo'' log
entries for uncommitted updates must be on disk before the page can be
stolen (written back).  On recovery, the LSN on the page reveals which
UNDO entries to apply to roll back the page. We use the absence of
commit records to figure out which transactions to roll back.

Thus, the single-page transactions of \yad work as follows.  An {\em
operation} consists of both a redo and an undo function, both of which
take one argument. An update is always the redo function applied to
the page (there is no ``do'' function), and it always ensures that the
redo log entry (with its LSN and argument) reaches the disk before
commit.  Similarly, an undo log entry, with its LSN and argument,
always reaches the disk before a page is stolen.  ARIES works
essentially the same way, but hard-codes recommended page 
formats and index structures~\cite{ariesIM}.

To manually abort a transaction, \yad could either reload the page
from disk and roll it forward to reflect committed transactions (this would imply ``no steal''), or it
could roll back the page using the undo entries applied in reverse LSN
order. (It currently does the latter.)


\eat{
Write-ahead-logging algorithms are quite simple if each operation
applied to the page file can be applied atomically.  This section will
describe a write ahead logging scheme that can transactionally update
a single page of storage that is guaranteed to be written to disk
atomically.  We refer the readers to the large body of literature
discussing write ahead logging if more detail is required.  Also, for
brevity, this section glosses over many standard write ahead logging
optimizations that \yad implements.


Assume an application wishes to transactionally apply a series of
functions to a piece of persistent storage.  For simplicity, we will
assume we have two deterministic functions, {\em undo}, and {\em
redo}.  Both functions take the contents of a page and a second
argument, and return a modified page.

As long as their second arguments match, undo and redo are inverses of
each other.  Normally, only calls to abort and recovery will invoke undo, so
we will assume that transactions consist of repeated applications of
the redo function.

Following the lead of ARIES (the write-ahead-logging system \yad
originally set out to implement), assume that the function is also
passed a distinct, monotonically increasing number each time it is
invoked, and that it records that number in an LSN (log sequence number)
field of the page.  In section~\ref{lsnFree}, we do away with this requirement.

We assume that while undo and redo are being executed, the
page they are modifying is pinned in memory.  Between invocations of
the two functions, the write-ahead-logging system may write the page
back to disk.  Also, multiple transactions may be interleaved, but
undo and redo must be executed atomically.  (However, \yad supports concurrent execution of operations.)

Finally, we assume that each invocation of redo and undo is recorded
in the log, along with a transaction id, LSN, and the argument passed into the redo or undo function.
(For efficiency, the page contents are not stored in the log.)

If abort is called during normal operation, the system will iterate
backwards over the log, invoking undo once for each invocation of redo
performed by the aborted transaction.  It should be clear that, in the
single transaction case, abort will restore the page to the state it
was in before the transaction began.  Note that each call to undo is
assigned a new LSN so the page LSN will be different.  Also, each undo
is also written to the log.
}

This section very briefly described how a simplified
write-ahead-logging algorithm might work, and glossed over many
details.  Like ARIES, \yad actually implements recovery in three
phases: Analysis, Redo and Undo.  

%Recovery is handled by playing the log forward, and only applying log
%entries that are newer than the version of the page on disk.  Once the
%end of the log is reached, recovery proceeds to abort any transactions
%that did not commit before the system crashed.\endnote{Like ARIES,
%\yad actually implements recovery in three phases, Analysis, Redo and
%Undo.}  Recovery arranges to continue any outstanding aborts where
%they left off, instead of rolling back the abort, only to restart it
%again.

\eat{
Note that recovery relies on the fact that it knows which version of
the page is recorded on disk, and that the page itself is
self-consistent.  If it passes an unknown version of a page into undo
(which is an arbitrary function), it has no way of predicting what
will happen.
}


\subsection{Multi-page transactions}

Of course, in practice, we wish to support transactions that span more
than one page.  Given a no-force/steal single-page transaction, this
is relatively easy.

First, we need to ensure that all log entries have a transaction ID
 so that we can tell that updates to different pages are part of
the same transaction (we need this in the single page case as well).
  Given single-page recovery, we can just apply it to
all of the pages touched by a transaction to recover a multi-page
transaction.  This works because steal and no-force already imply
that pages can be written back early or late (respectively), so there
is no need to write a group of pages back atomically.  In fact, we
need only ensure that redo entries for all pages reach the disk before
the commit record (and before commit returns).

\eat{
\subsection{Write-ahead-logging invariants}

In order to support recovery, a write-ahead-logging algorithm must
identify pages that {\em may} be written back to disk, and those that
{\em must} be written back to disk.  \yad provides full support for
Steal/no-Force write-ahead-logging, due to its generally favorable
performance properties.  ``Steal'' refers to the fact that pages may
be written back to disk before a transaction completes.  ``No-Force''
means that a transaction may commit before the pages it modified are
written back to disk.  

In a Steal/no-Force system, a page may be written to disk once the log
entries corresponding to the updates it contains are written to the
log file.  A page must be written to disk if the log file is full, and
the version of the page on disk is so old that deleting the beginning
of the log would lose redo information that may be needed at recovery.

Steal is desirable because it allows a single transaction to modify
more data than is present in memory.  Also, it provides more
opportunities for the buffer manager to write pages back to disk.
Otherwise, in the face of concurrent transactions that all modify the
same page, it may never be legal to write the page back to disk.  Of
course, if these problems would never come up in practice, an
application could opt for a no-Steal policy, possibly allowing it to
write less undo information to the log file.

No-Force is often desirable for two reasons.  First, forcing pages
modified by a transaction to disk can be extremely slow if the updates
are not near each other on disk.  Second, if many transactions update
a page, Force could cause that page to be written once for each transaction
that touched the page.  However, a Force policy could reduce the
amount of redo information that must be written to the log file.
}


\subsection{Nested top actions}
\label{sec:nta}
So far, we have glossed over the behavior of our system when concurrent
transactions modify the same data structure.  To understand the problems that
arise in this case, consider what
would happen if one transaction, A, rearranged the layout of a data
structure.  Next, assume a second transaction, B, modified that
structure, and then A aborted.  When A rolls back, its UNDO entries
will undo the rearrangement that it made to the data structure, without
regard to B's modifications.  This is likely to cause corruption.

Two common solutions to this problem are {\em total isolation} and
{\em nested top actions}.  Total isolation simply prevents any
transaction from accessing a data structure that has been modified by
another in-progress transaction.  An application can achieve this
using its own concurrency control mechanisms, or by holding a lock on
each data structure until the end of the transaction.  Releasing the
lock after the modification, but before the end of the transaction,
increases concurrency.  However, it means that follow-on transactions that use
that data may need to abort if a current transaction aborts ({\em
cascading aborts}).  %Related issues are studied in great detail in terms of optimistic concurrency control~\cite{optimisticConcurrencyControl, optimisticConcurrencyPerformance}.

Unfortunately, the long locks held by total isolation cause bottlenecks when applied to key
data structures.
Nested top actions are essentially mini-transactions that can
commit even if their containing transaction aborts; thus follow-on
transactions can use the data structure without fear of cascading
aborts.  

The key idea is to distinguish between the {\em logical operations} of a
data structure, such as inserting a key, and the {\em physical operations}
such as splitting tree nodes or or rebalancing a tree. The physical
operations do not need to be undone if the containing logical operation
(insert) aborts.  \diff{We record such operations using {\em logical 
logging} and {\em physical logging}, respectively.}

\diff{Each nested top action performs a single logical operation by applying
a number of physical operations to the page file.  Physical REDO log
entries are stored in the log so that recovery can repair any
temporary inconsistency that the nested top action introduces.
Logical UNDO entries are recorded so that the nested top action can be
rolled back even if concurrent transactions manipulate the data
structure.  Finally, physical UNDO entries are recorded so that 
the nested top action may be rolled back if the system crashes before 
it completes.}

\diff{When making use of nested top actions, we think of them as a
special type of latch that hides temporary inconsistencies from the
procedures executed during recovery.  Generally, such inconsistencies
must be hidden from other transactions in a multithreaded environment;
therefore we usually protect nested top actions with a mutex.}

\diff{This observation leads to the following mechanical conversion of
non-concurrent operations to thread-safe code that handles concurrent
transactions correctly:}

%Because nested top actions are easy to use and do not lead to 
%deadlock, we wrote a simple \yad extension that
%implements nested top actions.  The extension may be used as follows:

\begin{enumerate}
\item Wrap a mutex around each operation.  With care, it may be possible to use finer-grained locks, but it is rarely necessary.
\item Define a {\em logical} UNDO for each operation (rather than just
  using a set of page-level UNDO's).  For example, this is easy for a
  hashtable: the UNDO for {\em insert} is {\em remove}.  \diff{This logical
  undo function should arrange to acquire the mutex when invoked by
  abort or recovery.}
\item Add a ``begin nested
  top action'' right after the mutex acquisition, and a ``commit
  nested top action'' right before the mutex is released.  \diff{\yad provides a default nested top action implementation as an extension.}
\end{enumerate}

\noindent If the transaction that encloses the operation aborts, the logical
undo will {\em compensate} for its effects, leaving the structural
changes intact. 
% Note that this recipe does not ensure ISO transactional
%consistency and is largely orthogonal to the use of a lock manager.

We have found that it is easy to protect operations that make
structural changes to data structures with this recipe.
Therefore, we use them throughout our default data structure
implementations, although \yad does not preclude the use of more
complex schemes that lead to higher concurrency.


\subsection{Blind Writes}
\label{sec:blindWrites}
As described above, and in all database implementations of which we
are aware, transactional pages use LSNs on each page.  This makes it
difficult to map large objects onto multiple pages, as the LSNs break
up the object.  It is tempting to try to move the LSNs elsewhere, but
then they would not be written atomically with their page, which
defeats their purpose.

LSNs were introduced to prevent recovery from applying updates more
than once.  \diff{However, \yad can eliminate the LSN on each page by
constraining itself to deterministic REDO log entries that do not read
the contents of the page they update.}

%However, by constraining itself to a special type of idempotent redo and undo
%entries,\endnote{Idempotency does not guarantee that $f(g(x)) =
%  f(g(f(g(x))))$.  Therefore, idempotency does not guarantee that it is safe
%  to assume that a page is older than it is.}
%\yad can eliminate the LSN on each page.

Consider purely physical logging operations that overwrite a fixed
byte range on the page regardless of the page's initial state.  
We say that such operations perform ``blind writes.''
If all
operations that modify a page have this property, then we can remove
the LSN field, and have recovery \diff{use a conservative estimate 
of the LSN of each page that it is dealing with.}

\diff{For example, it 
could use the LSN of the most recent truncation point in the log, 
or during normal operation, \yad could occasionally write the 
LSN of the oldest dirty page to the log.}

% conservatively assume that it is
%dealing with a version of the page that is at least as old as the one
%on disk.  

To understand why this works, note that the log entries
update some subset of the bits on the page.  If the log entries do not
update a bit, then its value was correct before recovery began, so it
must be correct after recovery.  Otherwise, we know that recovery will
update the bit.  Furthermore, after all REDOs, the bit's value will be the
last value it contained before the crash, so we know that undo will behave
properly.

We call such pages ``LSN-free'' pages.  Although this technique is
novel for databases, it resembles the mechanism used by
RVM~\cite{lrvm}; \yad generalizes the concept and allows it to
co-exist with traditional pages.  Furthermore, efficient recovery and
log truncation require only minor modifications to our recovery
algorithm.  In practice, this is implemented by providing a buffer manager callback
for LSN free pages.  The callback computes a
conservative estimate of the page's LSN whenever the page is read from disk.
For a less conservative estimate, it suffices to write a page's LSN to
the log shortly after the page itself is written out; on recovery the
log entry is thus a conservative but close estimate.

Section~\ref{sec:zeroCopy} explains how LSN-free pages led us to new 
approaches for recoverable virtual memory and for large object storage.  
Section~\ref{sec:oasys} uses blind writes to efficiently update records 
on pages that are manipulated using more general operations.  \diff{We 
have not yet implemented LSN-free pages, so our experimental setup mimics
their behavior.}

\diff{Also note that while LSN-free pages assume that only bits that
are being updated will change, they do not assume that disk writes are
atomic.  Most disks do not atomically update more a single 512-byte
sector at a time.  However, most database systems make use of pages
that are larger than 512 bytes.  Recovery schemes that rely upon LSN
fields in pages must detect and deal with torn pages
directly~\cite{tornPageStuffMohan}.  Because LSN-free page recovery
does not assume page writes are atomic, it handles torn pages with no
extra effort.}


\subsection{Media recovery}

\diff{Hard drives may lose data due to hardware failures, or because a
sector is being written when power is lost.  The drive hardware stores a
checksum with each sector, and will issue a read error if the checksum
does not match~\cite{something}.}  Like ARIES, \yad can recover lost pages in the page
file by reinitializing the page to zero, and playing back the entire
log.  In practice, a system administrator would periodically back up
the page file up, thus enabling log truncation and shortening recovery
time.

\eat{  This is pretty redundant.
\subsection{Modular operations semantics}

The smallest unit of a \yad transaction is the {\em operation}.  An
operation consists of a {\em redo} function, {\em undo} function, and
a log format.  At runtime or if recovery decides to reapply the
operation, the redo function is invoked with the contents of the log
entry as an argument.  During abort, or if recovery decides to undo
the operation, the undo function is invoked with the contents of the
log as an argument.  Like Berkeley DB, and most database toolkits, we
allow system designers to define new operations.  Unlike earlier
systems, we have based our library of operations on object oriented
collection libraries, and have built complex index structures from
simpler structures.  These modules are all directly available,
providing a wide range of data structures to applications, and
facilitating the develop of more complex structures through reuse.  We
compare the performance of our modular approach with a monolithic
implementation on top of \yad, using Berkeley DB as a baseline.
}

\eat{ \subsection{Buffer manager policy}

Generally, write ahead logging algorithms ensure that the most recent
version of each memory-resident page is stored in the buffer manager,
and the most recent version of other pages is stored in the page file.
This allows the buffer manager to present a uniform view of the stored
data to the application.  The buffer manager uses a cache replacement
policy (\yad currently uses LRU-2 by default) to decide which pages
should be written back to disk.

Section~\ref{sec:oasys}, we will provide example where the most recent
version of application data is not managed by \yad at all, and
Section~\ref{zeroCopy} explains why efficiency may force certain
operations to bypass the buffer manager entirely.


\subsection{Durability}

\eat{\yad makes use of the same basic recovery strategy as existing
write-ahead-logging schemes such as ARIES.  Recovery consists of three
stages, {\em analysis}, {\em redo}, and {\em undo}.  Analysis is
essentially a performance optimization, and makes use of information
left during forward operation to reduce the cost of redo and undo.  It
also decides which transactions committed, and which aborted.  The
redo phase iterates over the log, applying the redo function of each
logged operation if necessary.  Once the log has been played forward,
the page file and buffer manager are in the same conceptual state they
were in at crash.  The undo phase simply aborts each transaction that
does not have a commit entry, exactly as it would during normal
operation.
}
%From the application's perspective, logging and durability are interesting for a
%number of reasons.  First, 
If full transactional durability is
unneeded, the log can be flushed to disk less frequently, improving
performance.  In fact, \yad allows applications to store the
transaction log in memory, reducing disk activity at the expense of
recovery.  We are in the process of optimizing the system to handle
fully in-memory workloads efficiently.  Of course, durability is closely
tied to system management issues such as reliability, replication and so on.  
These issues are beyond the scope of this discussion.  Section~\ref{logReordering} will describe why applications might decide to manipulate the log directly.
}
\subsection{Summary of Transactional Pages}

This section provided an extremely brief overview of transactional
pages and write-ahead-logging.  Transactional pages are a valuable
building block for a wide variety of data management systems, as we
show in the next section.  Nested top actions and LSN-free pages
enable important optimizations.  In particular, \yad allows general 
custom operations using LSNs, or custom blind-write operations
without LSNs.  This enables transactional manipulation of large, 
contiguously stored objects.

\eat{
Although the extensions that it proposes
require a fair amount of knowledge about transactional logging
schemes, our initial experience customizing the system for various
applications is positive.  We believe that the time spent customizing
the library is less than amount of time that it would take to work
around typical problems with existing transactional storage systems.

%However, we do not yet have a good understanding of the practical testing and
%reliability issues that arise as the system is modified in
%this fashion.
}



\section{Extending \yad}
\label{sec:extensions}

\diff{The previous section described how \yad implements conventional
transactional storage.  In this section we discuss ways in which \yad
can be customized to provide more specialized transactions.  First,
the mechanisms that allow new operations will be discussed.  These
mechanisms provide the base of \yads customizable page formats and
ability to support application-specific transactional data structures.
Next, an example of how \yads recovery mechanism can be changed will
be discussed.}

\diff{In this section we break some of the typical assumptions made by
transactional storage algorithms.  The discussion of custom log
operations updates pages at the byte level, and describes how one
might implement functions that organize pages into records, or to
provide more exotic semantics.}

\diff{The customized recovery algorithm removes LSN's from pages, and
instead opts to estimate LSN's during recovery, and recalculate them
during normal forward operation.  This in turn breaks the reliance on
pages as an atomic unit of recovery, but prevents us from using most
conventional database page layout techniques.}

\diff{This section discusses changes that are made at multiple levels
of abstraction.  We will attempt to describe which level is being
described, and the semantics provided by the levels it builds upon.}

%This section describes proof-of-concept extensions to \yad.
%Performance figures accompany the extensions that we have implemented.
%We discuss existing approaches to the systems presented here when
%appropriate.

\subsection{Adding log operations}
\rcs{This section needs to be merged into the new section 3, because that is where we discuss how to add new log operations.  (In with the new nested top action stuff, probably).  That will leave a section to focus on LSN-free pages, and other things that break the ARIES assumptions.  That way, blind writes and lsn-free pages can be in the same place.}
\label{sec:wal}
\begin{figure}
\includegraphics[%
   width=1\columnwidth]{figs/structure.pdf}
\caption{\sf\label{fig:structure} The portions of \yad that directly interact with new operations.}
\end{figure}
\yad allows application developers to easily add new operations to the
system.  Many of the customizations described below can be implemented
using custom log operations.  In this section, we describe how to implement an
``ARIES style'' concurrent, steal/no-force operation using 
\diff{physical redo, logical undo} and per-page LSN's.
Such operations are typical of high-performance commercial database
engines.

As we mentioned above, \yad operations must implement a number of
functions.  Figure~\ref{fig:structure} describes the environment that
schedules and invokes these functions.  The first step in implementing
a new set of log interfaces is to decide upon an interface that these log
interfaces will export to callers outside of \yad.  

The externally visible interface is implemented by wrapper functions
and read-only access methods.  The wrapper function modifies the state
of the page file by packaging the information that will be needed for
undo and redo into a data format of its choosing.  This data structure
is passed into Tupdate().  Tupdate() copies the data to the log, and
then passes the data into the operation's REDO function.
 
REDO modifies the page file directly (or takes some other action).  It
is essentially an interpreter for the log entries it is associated
with.  UNDO works analogously, but is invoked when an operation must
be undone (usually due to an aborted transaction, or during recovery).

This pattern applies in many cases.  In
order to implement a ``typical'' operation, the operations
implementation must obey a few more invariants:

\begin{itemize}
\item Pages should only be updated inside REDO and UNDO functions.
\item Page updates atomically update the page's LSN by pinning the page.
\item If the data seen by a wrapper function must match data seen
  during REDO, then the wrapper should use a latch to protect against
  concurrent attempts to update the sensitive data (and against
  concurrent attempts to allocate log entries that update the data).
\item Nested top actions (and logical undo), or ``big locks'' (total isolation but lower concurrency) should be used to implement multi-page updates. (Section~\ref{sec:nta})
\end{itemize}

\subsection{LSN-Free pages}
\label{sec:zeroCopy}
In Section~\ref{sec:blindWrites}, we describe how operations can avoid recording
LSN's on the pages they modify.  Essentially, operations that update pages \diff{without examining their contents}
% make use of purely physical logging 
need not heed page boundaries.
%, as physiological operations must.  
Recall that purely physical logging
interacts poorly with concurrent transactions that modify the same
data structures or pages, so LSN-Free pages are not applicable in all
situations. \rcs{I think we can support physiological logging; once REDO is done, we know the LSN.  Why not do logical UNDO?}

Consider the retrieval of a large (page spanning) object stored on
pages that contain LSN's.  The object's data will not be contiguous.
Therefore, in order to retrieve the object, the transaction system must
load the pages contained on disk into memory, and perform a byte-by-byte copy of the
portions of the pages that contain the large object's data into a second buffer.  

Compare
this approach to a modern filesystem, which allows applications to
perform a DMA copy of the data into memory, avoiding the expensive
byte-by-byte copy of the data, and allowing the CPU to be used for
more productive purposes.  Furthermore, modern operating systems allow
network services to use DMA and network adaptor hardware to read data
from disk, and send it over a network socket without passing it
through the CPU.  Again, this frees the CPU, allowing it to perform
other tasks.

We believe that LSN free pages will allow reads to make use of such
optimizations in a straightforward fashion.  Zero copy writes are more challenging, but could be
performed by performing a DMA write to a portion of the log file.
However, doing this complicates log truncation, and does not address
the problem of updating the page file.  We suspect that contributions
from the log based filesystem~\cite{lfs} literature can address these problems in
a straightforward fashion.  In particular, we imagine storing 
portions of the log (the portion that stores the blob) in the 
page file, or other addressable storage.  In the worst case, 
the blob would have to be relocated in order to defragment the 
storage.  Assuming the blob was relocated once, this would amount 
to a total of three, mostly sequential disk operations.  (Two 
writes and one read.)  However, in the best case, the blob would only need to written once.
In contrast, a conventional atomic blob implementation would always need 
to write the blob twice. %but also may need to create complex 
%structures such as B-Trees, or may evict a large number of 
%unrelated pages from the buffer pool as the blob is being written 
%to disk.  

Alternatively, we could use DMA to overwrite the blob in the page file
in a non-atomic fashion, providing filesystem style semantics.
(Existing database servers often provide this mode based on the
observation that many blobs are static data that does not really need
to be updated transactionally.~\cite{sqlserver}) Of course, \yad could
also support other approaches to blob storage, such as B-Tree layouts
that allow arbitrary insertions and deletions in the middle of
objects~\cite{esm}.

Finally, RVM, recoverable virtual memory, made use of LSN-free pages
so that it could use mmap() to map portions of the page file into
application memory\cite{lrvm}.  However, without support for logical log entries
and nested top actions, it would be difficult to implement a
concurrent, durable data structure using RVM.  We plan to add RVM
style transactional memory to \yad in a way that is compatible with
fully concurrent collections such as hash tables and tree structures.

\section{Experiments}
\subsection{Experimental setup}



\label{sec:experimental_setup}

We chose Berkeley DB in the following experiments because, among
commonly used systems, it provides transactional storage primitives
that are most similar to \yad.  Also, Berkeley DB is designed to provide high
performance and high concurrency.  For all tests, the two libraries
provide the same transactional semantics, unless explicitly noted.

All benchmarks were run on an Intel Xeon 2.8 GHz with 1GB of RAM and a
10K RPM SCSI drive formatted using with ReiserFS~\cite{reiserfs}.\endnote{We found that the
  relative performance of Berkeley DB and \yad under single threaded testing is sensitive to
  filesystem choice, and we plan to investigate the reasons why the
  performance of \yad under ext3 is degraded. However, the results
  relating to the \yad optimizations are consistent across filesystem
  types.}  All results correspond to the mean of multiple runs with a
95\% confidence interval with a half-width of 5\%.

We used Berkeley DB 4.2.52 as it existed in Debian Linux's testing
branch during March of 2005, with the flags DB\_TXN\_SYNC, and
DB\_THREAD enabled. These flags were chosen to match Berkeley DB's
configuration to \yads as closely as possible.  In cases where
Berkeley DB implements a feature that is not provided by \yad, we
only enable the feature if it improves Berkeley DB's performance.

Optimizations to Berkeley DB that we performed included disabling the
lock manager, though we still use ``Free Threaded'' handles for all
tests.  This yielded a significant increase in performance because it
removed the possibility of transaction deadlock, abort, and
repetition.  However, disabling the lock manager caused highly
concurrent Berkeley DB benchmarks to become unstable, suggesting either a
bug or misuse of the feature.  

With the lock manager enabled, Berkeley
DB's performance in the multithreaded test in Section~\ref{sec:lht} strictly decreased with
increased concurrency.  (The other tests were single-threaded.)  We also
increased Berkeley DB's buffer cache and log buffer sizes to match
\yads default sizes.

We expended a considerable effort tuning Berkeley DB, and our efforts
significantly improved Berkeley DB's performance on these tests.
Although further tuning by Berkeley DB experts would probably improve
Berkeley DB's numbers, we think that we have produced a reasonably
fair comparison.  The results presented here have been reproduced on
multiple machines and file systems.

\subsection{Linear hash table}
\label{sec:lht}
\begin{figure}[t]
\includegraphics[%
   width=1\columnwidth]{figs/bulk-load.pdf}
%\includegraphics[%
%   width=1\columnwidth]{bulk-load-raw.pdf}
%\vspace{-30pt}
\caption{\sf\label{fig:BULK_LOAD} Performance of \yad and Berkeley DB hashtable implementations.  The
test is run as a single transaction, minimizing overheads due to synchronous log writes.}
\end{figure}
\begin{figure}[t]
%\hspace*{18pt}
%\includegraphics[%
%   width=1\columnwidth]{tps-new.pdf}
\includegraphics[%
   width=1\columnwidth]{figs/tps-extended.pdf}
%\vspace{-36pt}
\caption{\sf\label{fig:TPS} High concurrency performance of Berkeley DB and \yad.  We were unable to get Berkeley DB to work correctly with more than 50 threads. (See text)
}
\end{figure}

Although the beginning of this paper describes the limitations of
physical database models and relational storage systems in great
detail, these systems are the basis of most common transactional
storage routines.  Therefore, we implement a key-based access 
method in this section. We argue that 
obtaining reasonable performance in such a system under \yad is
straightforward.  We then compare our simple, straightforward 
implementation to our hand-tuned version and Berkeley DB's implementation.

The simple hash table uses nested top actions to atomically update its
internal structure.  It uses a {\em linear} hash function~\cite{lht}, allowing
it to incrementally grow its buffer list.  It is based on a number of
modular subcomponents.  Notably, its bucket list is a growable array
of fixed length entries (a linkset, in the terms of the physical
database model) and the user's choice of two different linked list
implementations.

The hand-tuned hashtable also uses a linear hash
function.  However, it is monolithic and uses carefully ordered writes to
reduce runtime overheads such as log bandwidth.  Berkeley DB's
hashtable is a popular, commonly deployed implementation, and serves
as a baseline for our experiments.

Both of our hashtables outperform Berkeley DB on a workload that
bulk loads the tables by repeatedly inserting (key, value) pairs.
%although we do not wish to imply this is always the case.
%We do not claim that our partial implementation of \yad
%generally outperforms, or is a robust alternative
%to Berkeley DB.  Instead, this test shows that \yad is comparable to
%existing systems, and that its modular design does not introduce gross
%inefficiencies at runtime.
The comparison between the \yad  implementations is more
enlightening.  The performance of the simple hash table shows that
straightforward data structure implementations composed from
simpler structures can perform as well as the implementations included 
in existing monolithic systems.  The hand-tuned
implementation shows that \yad allows application developers to
optimize key primitives.

% I cut this because Berkeley db supports custom data structures....

%In the
%best case, past systems allowed application developers to provide
%hints to improve performance.  In the worst case, a developer would be
%forced to redesign and application to avoid sub-optimal properties of
%the transactional data structure implementation.

Figure~\ref{fig:TPS} describes the performance of the two systems under
highly concurrent workloads.  For this test, we used the simple
(unoptimized) hash table, since we are interested in the performance of a
clean, modular data structure that a typical system implementor might
 produce, not the performance of our own highly tuned,
monolithic implementations.

Both Berkeley DB and \yad can service concurrent calls to commit with
a single synchronous I/O.\endnote{The multi-threaded benchmarks
  presented here were performed using an ext3 filesystem, as high
  concurrency caused both Berkeley DB and \yad to behave unpredictably
  when ReiserFS was used.  However, \yads multi-threaded throughput
  was significantly better that Berkeley DB's under both filesystems.}
\yad scaled quite well, delivering over 6000 transactions per
second,\endnote{The concurrency test was run without lock managers, and the
  transactions obeyed the A, C, and D properties.  Since each
  transaction performed exactly one hashtable write and no reads, they also
  obeyed I (isolation) in a trivial sense.}  and provided roughly
double Berkeley DB's throughput (up to 50 threads).  We do not report
the data here, but we implemented a simple load generator that makes
use of a fixed pool of threads with a fixed think time.  We found that
the latency of Berkeley DB and \yad were similar, showing that \yad is 
not simply trading latency for throughput during the concurrency benchmark.


\begin{figure*}
\includegraphics[width=1\columnwidth]{figs/object-diff.pdf}
\hspace{.2in}
\includegraphics[width=1\columnwidth]{figs/mem-pressure.pdf}
\vspace{-.15in}
\caption{\sf \label{fig:OASYS}
The effect of \yad object serialization optimizations under low and high memory pressure.}
\end{figure*}

\subsection{Object persistence}
\label{sec:oasys}
Numerous schemes are used for object serialization.  Support for two
different styles of object serialization have been implemented in
\yad.  We could have just as easily implemented a persistence
mechanism for a statically typed functional programming language, a
dynamically typed scripting language, or a particular application,
such as an email server.  In each case, \yads lack of a hard-coded data
model would allow us to choose the representation and transactional
semantics that make the most sense for the system at hand.

The first object persistence mechanism, pobj, provides transactional updates to objects in
Titanium, a Java variant.  It transparently loads and persists
entire graphs of objects, but will not be discussed in further detail.

The second variant was built on top of a C++ object
serialization library, \oasys.  \oasys makes use of pluggable storage
modules that implement persistent storage, and includes plugins
for Berkeley DB and MySQL.  

This section will describe how the \yad
\oasys plugin reduces amount of data written to log, while using half as much system
memory as the other two systems.

We present three variants of the \yad plugin here.  The first treats \yad like
Berkeley DB.  The second, ``update/flush'' customizes the behavior of the buffer
manager.  Instead of maintaining an up-to-date version of each object
in the buffer manager or page file, it allows the buffer manager's
view of live application objects to become stale.  This is safe since
the system is always able to reconstruct the appropriate page entry
from the live copy of the object.

By allowing the buffer manager to contain stale data, we reduce the
number of times the \yad \oasys plugin must update serialized objects in the buffer manager.
%  Reducing the number of serializations decreases
%CPU utilization, and it also
This allows us to drastically decrease the
size of the page file.  In turn this allows us to increase the size of
the application's cache of live objects.

We implemented the \yad buffer-pool optimization by adding two new
operations, update(), which only updates the log, and flush(), which
updates the page file.  

The reason it would be difficult to do this with Berkeley DB is that
we still need to generate log entries as the object is being updated.
Otherwise, commit would not be durable, unless we queued up log 
entries, and wrote them all before committing.  
  This would cause Berkeley DB to write data back to the
page file, increasing the working set of the program, and increasing
disk activity.

Furthermore, objects may be written to disk in an
order that differs from the order in which they were updated, 
violating one of the write-ahead-logging invariants.  One way to 
deal with this is to maintain multiple LSN's per page.  This means we would need to register a
callback with the recovery routine to process the LSN's (a similar
callback will be needed in Section~\ref{sec:zeroCopy}), and 
extend \yads page format to contain per-record LSN's.  
Also, we must prevent \yads storage allocation routine from overwriting the per-object 
LSN's of deleted objects that may still be addressed during abort or recovery.  

Alternatively, we could arrange for the object pool to cooperate 
further with the buffer pool by atomically updating the buffer 
manager's copy of all objects that share a given page, removing the 
need for multiple LSN's per page, and simplifying storage allocation.

However, the simplest solution, and the one we take here, is based on the observation that
updates (not allocations or deletions) of fixed length objects are blind writes.
This allows us to do away with per-object LSN's entirely.  Allocation and deletion can then be handled
as updates to normal LSN containing pages.  At recovery time, object
updates are executed based on the existence of the object on the page
and a conservative estimate of its LSN.  (If the page doesn't contain
the object during REDO then it must have been written back to disk
after the object was deleted.  Therefore, we do not need to apply the
REDO.)  This means that the system can ``forget'' about objects that 
were freed by committed transactions, simplifying space reuse 
tremendously.

The third \yad plugin, ``delta'' incorporates the buffer
manager optimizations.  However, it only writes the changed portions of
objects to the log.  Because of \yads support for custom log entry
formats, this optimization is straightforward.

%In addition to the buffer-pool optimizations, \yad provides several 
%options to handle UNDO records in the context
%of object serialization. The first is to use a single transaction for
%each object modification, avoiding the cost of generating or logging
%any UNDO records. The second option is to assume that the
%application will provide a custom UNDO for the delta, 
%which increases the size of the log entry generated by each update, 
%but still avoids the need to read or update the page
%file.
%
%The third option is to relax the atomicity requirements for a set of
%object updates and again avoid generating any UNDO records. This
%assumes that the application cannot abort individual updates, 
%and is willing to
%accept that some prefix of logged but uncommitted updates may 
%be applied to the page
%file after recovery. 

\oasys does not export transactions to its callers.  Instead, it
is designed to be used in systems that stream objects over an
unreliable network connection.  Each object update corresponds to an
independent message, so there is never any reason to roll back an
applied object update.  On the other hand, \oasys does support a
flush method, which guarantees the durability of updates after it
returns.  In order to match these semantics as closely as possible,
\yads update/flush and delta optimizations do not write any 
undo information to the log.  

These ``transactions'' are still durable
after commit, as commit forces the log to disk. 
%For the benchmarks below, we
%use this approach, as it is the most aggressive and is
As far as we can tell, MySQL and Berkeley DB do not support this
optimization in a straightforward fashion.  (``Auto-commit'' comes
close, but does not quite provide the correct durability semantics.)
%not supported by any other general-purpose transactional 
%storage system (that we know of).

The operations required for these two optimizations required 
150 lines of C code, including whitespace, comments and boilerplate
function registrations.\endnote{These figures do not include the
  simple LSN free object logic required for recovery, as \yad does not
  yet support LSN free operations.}  Although the reasoning required
to ensure the correctness of this code is complex, the simplicity of
the implementation is encouraging.

In this experiment, Berkeley DB was configured as described above.  We
ran MySQL using InnoDB for the table engine.  For this benchmark, it
is the fastest engine that provides similar durability to \yad. We
linked the benchmark's executable to the libmysqld daemon library,
bypassing the RPC layer. In experiments that used the RPC layer, test
completion times were orders of magnitude slower.

Figure~\ref{fig:OASYS} presents the performance of the three
\yad optimizations, and the \oasys plugins implemented on top of other
systems.  As we can see, \yad performs better than the baseline
systems, which is not surprising, since it is not providing the A
property of ACID transactions.  (Although it is applying each individual operation atomically.)

In non-memory bound systems, the optimizations nearly double \yads
performance by reducing the CPU overhead of object serialization and
the number of log entries written to disk.  In the memory bound test,
we see that update/flush indeed improves memory utilization.


\subsection{Manipulation of logical log entries}
\label{sec:logging}
\begin{figure}
\includegraphics[width=1\columnwidth]{figs/graph-traversal.pdf}
\vspace{-24pt}
\caption{\sf\label{fig:multiplexor} Because pages are independent, we
can reorder requests among different pages. Using a log demultiplexer,
we partition requests into independent queues, which can be 
handled in any order, improving locality and merging opportunities.}
\end{figure}
\begin{figure}[t]
\includegraphics[width=1\columnwidth]{figs/oo7.pdf}
\vspace{-15pt}
\caption{\sf\label{fig:oo7} oo7 benchmark style graph traversal.  The optimization performs well due to the presence of non-local nodes.}
\end{figure}

\begin{figure}[t]
\includegraphics[width=1\columnwidth]{figs/trans-closure-hotset.pdf}
\vspace{-12pt}
\caption{\sf\label{fig:hotGraph} Hot set based graph traversal for random graphs with out-degrees of 3 and 9.  Here
we see that the multiplexer helps when the graph has poor locality.
In the cases where depth first search performs well, the
reordering is inexpensive.}
\end{figure}

Database optimizers operate over relational algebra expressions that
correspond to logical operations over streams of data.  \yad
does not provide query languages, relational algebra, or other such query processing primitives.  

However, it does include an extensible logging infrastructure.
Furthermore, \diff{most operations that support concurrent transactions already
provide logical UNDO (and therefore logical REDO, if each operation has an
inverse).}
%many
%operations that make use of physiological logging implicitly
%implement UNDO (and often REDO) functions that interpret logical
%requests.

Logical operations often have some nice properties that this section
will exploit.  Because they can be invoked at arbitrary times in the
future, they tend to be independent of the database's physical state.
Often, they correspond to operations that programmers understand.

Because of this, application developers can easily determine whether
logical operations may be reordered, transformed, or even
dropped from the stream of requests that \yad is processing.

If requests can be partitioned in a natural way, load
balancing can be implemented by splitting requests across many nodes.
Similarly, a node can easily service streams of requests from multiple
nodes by combining them into a single log, and processing the log
using operation implementations.  For example, this type of optimization 
is used by RVM's log-merging operations~\cite{lrvm}.

Furthermore, application-specific
procedures that are analogous to standard relational algebra methods
(join, project and select) could be used to efficiently transform the data
while it is still layed out sequentially
in non-transactional memory.

%Note that read-only operations do not necessarily generate log
%entries.  Therefore, applications may need to implement custom
%operations to make use of the ideas in this section.

%Although \yad has rudimentary support for a \diff{cluster hash table\cite{cht}} that uses 
%two-phase commit to recover from node crashes}, we have not yet implemented networking primitives for logical logs.
\rcs{Cut sentence about two-phase commit cluster hash table, networking primitves for logical logs.}
Therefore, we implemented a single node log-reordering scheme that increases request locality
during the traversal of a random graph.  The graph traversal system
takes a sequence of (read) requests, and partitions them using some
function.  It then processes each partition in isolation from the
others.  We considered two partitioning functions.  The first divides the page file
into equally sized contiguous regions, which increases locality.  The second takes the hash
of the page's offset in the file, which enables load balancing.
%%  The second policy is interesting
%The first, partitions the
%requests according to the hash of the node id they refer to, and would be useful for load balancing over a network.
%(We expect the early phases of such a traversal to be bandwidth, not
%latency limited, as each node would stream large sequences of
%asynchronous requests to the other nodes.) 

Our benchmarks partition requests by location.  We chose the
position size so that each partition can fit in \yads buffer pool.

We ran two experiments.  Both stored a graph of fixed size objects in
the growable array implementation that is used as our linear
hashtable's bucket list.
The first experiment (Figure~\ref{fig:oo7})
is loosely based on the oo7 database benchmark.~\cite{oo7}.  We
hard-code the out-degree of each node, and use a directed graph.  OO7
constructs graphs by first connecting nodes together into a ring.
It then randomly adds edges between the nodes until the desired
out-degree is obtained.  This structure ensures graph connectivity.
If the nodes are laid out in ring order on disk then it also ensures that
one edge from each node has good locality while the others generally
have poor locality.

The second experiment explicitly measures the effect of graph locality
on our optimization (Figure~\ref{fig:hotGraph}). It extends the idea
of a hot set to graph generation.  Each node has a distinct hot set
that includes the 10\% of the nodes that are closest to it in ring
order.  The remaining nodes are in the cold set.  We use random edges
instead of ring edges for this test.  This does not ensure graph
connectivity, but we used the same random seeds for the two systems.

When the graph has good locality, a normal depth first search
traversal and the prioritized traversal both perform well.  The
prioritized traversal is slightly slower due to the overhead of extra
log manipulation. As locality decreases, the partitioned traversal
algorithm's outperforms the naive traversal.


\section{Related Work}

This paper has described a number of custom transactional storage
extensions, and explained why can \yad support them.  This section
will describe existing ideas in the literature that we would like to
incorporate into \yad.  An overview of database systems that have 
goals similar to our own is in Section~\ref{sec:otherDBs}.

Different large object storage systems provide different API's.
Some allow arbitrary insertion and deletion of bytes~\cite{esm} or
pages~\cite{sqlserver} within the object, while typical filesystems
provide append-only storage allocation~\cite{ffs}.
Record-oriented file systems are an older, but still-used~\cite{gfs}
alternative. Each of these API's addresses 
different workloads.

Although most filesystems attempt to lay out data in logically sequential
order, write-optimized filesystems lay files out in the order they
were written~\cite{lfs}.  Schemes to improve locality between small
objects exist as well. Relational databases allow users to specify the order
in which tuples will be layed out, and often leave portions of pages
unallocated to reduce fragmentation as new records are allocated.

Memory allocation routines also address this problem.  For example, the Hoard memory
allocator is a highly concurrent version of malloc that
makes use of thread context to allocate memory in a way that favors
cache locality~\cite{hoard}.  %Other work makes use of the caller's stack to infer
%information about memory management.~\cite{xxx} \rcs{Eric, do you have
%  a reference for this?}

Finally, many systems take a hybrid approach to allocation.  Examples include
databases with blob support, and a number of
filesystems~\cite{reiserfs,ffs}.

We are interested in allowing applications to store records in
the transaction log.  Assuming log fragmentation is kept to a
minimum, this is particularly attractive on a single disk system.  We
plan to use ideas from LFS~\cite{lfs} and POSTGRES~\cite{postgres}
to implement this.

Starburst~\cite{starburst} provides a flexible approach to index
management, and database trigger support, as well as hints for small
object layout.

The Boxwood system provides a networked, fault-tolerant transactional
B-Tree and ``Chunk Manager.''  We believe that \yad is an interesting
complement to such a system, especially given \yads focus on
intelligence and optimizations within a single node, and Boxwood's
focus on multiple node systems.  In particular, it would be
interesting to explore extensions to the Boxwood approach that make
use of \yads customizable semantics (Section~\ref{sec:wal}), and fully logical logging
mechanism. (Section~\ref{sec:logging})

\section{Future Work}

Complexity problems may begin to arise as we attempt to implement more
extensions to \yad.  However, \yads implementation is still fairly simple:

\begin{itemize}
\item The core of \yad is roughly 3000 lines
of C code, and implements the buffer manager, IO, recovery, and other
systems
\item Custom operations account for another 3000 lines of code
\item Page layouts and logging implementations account for 1600 lines of code.
\end{itemize}

The complexity of the core of \yad is our primary concern, as it
contains the hard-coded policies and assumptions.  Over time, the core has
shrunk as functionality has been moved into extensions.  We expect
this trend to continue as development progresses.  

A resource manager
is a common pattern in system software design, and manages
dependencies and ordering constraints between sets of components.
Over time, we hope to shrink \yads core to the point where it is
simply a resource manager and a set of implementations of a few unavoidable
algorithms related to write-ahead-logging.  For instance, 
we suspect that support for appropriate callbacks will 
allow us to hard-code a generic recovery algorithm into the 
system.  Similarly, any code that manages book-keeping information, such as 
LSN's may be general enough to be hard-coded.  

Of course, we also plan to provide \yads current functionality, including the algorithms
mentioned above as modular, well-tested extensions.
Highly specialized \yad extensions, and other systems would be built
by reusing \yads default extensions and implementing new ones.


\section{Conclusion}

We have presented \yad, a transactional storage library that addresses
the needs of system developers.  \yad provides more opportunities for
specialization than existing systems.  The effort required to extend
\yad to support a new type of system is reasonable, especially when
compared to currently common practices, such as working around
limitations of existing systems, breaking guarantees regarding data
integrity, or reimplementing the entire storage infrastructure from
scratch.

We have demonstrated that \yad provides fully
concurrent, high performance transactions, and explained how it can
support a number of systems that currently make use of suboptimal or
ad-hoc storage approaches.  Finally, we have explained how \yad can be
extended in the future to support a larger range of systems.

\section{Acknowledgements}

The idea behind the \oasys buffer manager optimization is from Mike
Demmer.  He and Bowei Du implemented \oasys.  Gilad Arnold and Amir Kamil implemented
 for pobj.  Jim Blomo, Jason Bayer, and Jimmy
Kittiyachavalit worked on an early version of \yad.

Thanks to C. Mohan for pointing out the need for tombstones with
per-object LSN's.  Jim Gray provided feedback on an earlier version of
this paper, and suggested we use a resource manager to manage
dependencies within \yads API.  Joe Hellerstein and Mike Franklin
provided us with invaluable feedback.

\section{Availability}

Additional information, and \yads source code is available at:

\begin{center}
%{\tt http://www.cs.berkeley.edu/sears/\yad/}
{\small{\tt http://www.cs.berkeley.edu/\ensuremath{\sim}sears/\yad/}}
%{\tt http://www.cs.berkeley.edu/sears/\yad/}
\end{center}

{\footnotesize \bibliographystyle{acm}
\nocite{*}
\bibliography{LLADD}}

\theendnotes

\end{document}







